Public key cryptography using matrices

ABSTRACT

The invention provides techniques for secure messages transmission using a public key system to exchange secret keys. A first entity creates public and private keys by generating a product n of two large, randomly chosen prime numbers, and then generating random matrices {A, C}, in the group GL(r,Z n ) with a chosen matrix rank r such that AC is not equal to CA, and then generating a matrix B=CAC, and finding a matrix G that commutes with C. Matrices A, B, G and the integers n and r are then published as the public key and matrix C is then kept as the private key. A second entity then obtains the public key and calculates a secret matrix D that commutes with G, and further calculates the matrices K=DBD and E=DAD. The message to be sent is then encrypted using matrix K as the secret key and then sent to the first entity with matrix E. First entity then retrieves secret matrix K using K=CEC and then decrypts the received encrypted message using the retrieved secret matrix K.

RELATED APPLICATIONS

This application is a Divisional of U.S. application Ser. No.10/260,818, filed Sep. 30, 2002 now U.S. Pat. No. 7,184,551, which isincorporated herein by reference.

FIELD OF THE INVENTION

The present invention relates to data communications, and moreparticularly to cryptography for secure data transmission.

BACKGROUND OF THE INVENTION

Electronic messages are generally transmitted between remotecorrespondents via a communications system typically including a networkof interconnected computers. Such messages are readily intercepted andviewed by others using the network. Thus, correspondents desiringprivacy may encrypt or encode a message such that only the recipient candecrypt or decode the message to view the contents.

In a public key encryption system, a person wishing to receive encryptedmessages (a potential recipient) is able to generate a special set ofnumeric values. Some of these numeric values are published by therecipient as a public key and the remaining numeric values are kept asthe recipient's private key. A second person (a sender) wishing to sendan encrypted message to the recipient, first obtains the recipient'spublic key, and then encrypts a message using this public keyinformation. The message is then sent to the recipient. The recipient isthen able to use his or her private key information to decrypt theencrypted message much more rapidly than a message eavesdropper who doesnot have the private key information. In all public key schemes known,there is a mathematical relationship between the private key and thepublic key. Finding the private key via the mathematical relationshipcan be made arbitrarily difficult at the expense of encryption and/ordecryption performance.

A well-known encryption technique is disclosed in U.S. Pat. No.4,405,829 to Rivest et al., which is incorporated by reference. Thetechnique is also known as the RSA public key system. The RSA algorithmperforms integer arithmetic modulo n, where n is a product of two large,randomly chosen prime numbers. A recipient generates a private exponentkey using knowledge of the prime factors and a chosen public exponent.The public exponent and modulus n is published as the public key. Themessage sender uses the public key information to break up messages intopieces, each of which is numerically encoded in an agreed-on format tolie in the modulus range. The sender then takes each piece of themessage as a numeric value and raises it to the public exponent, withthe result calculated as modulo n. The result of encoding each piece isan encrypted value.

The above-described “power-mod” process is generally fast for smallpowers, so public exponents, tend to be relatively small compared to n.The sender then packs all the values in an agreed-on format to form theencrypted message. The recipient takes the message and breaks it up intothe same sets of encrypted values modulo n. For each value, therecipient raises the encrypted message to their private exponent modulon. This results in using the power-mod function again. Each resultingvalue is then unpacked to reclaim the original encrypted message.

To ensure security, n must be chosen so that factorization into itsprime factors is not feasible using the fastest known algorithms. If n'sfactors can be found, then the private exponent can be easilycalculated. Unfortunately, in terms of performance, the private exponentis generally a large number less than the modulus n, and the power-modfunction is relatively slow for large n when compared withmultiplication.

For a secure 1024-bit modulus n, a typical 1 GHz processor can encryptdata using the RSA algorithm with a secure public exponent of 2¹⁶+1 at arate of around 125,000 bits per second. Decryption is around 50 timesslower at about 2,500 bits per second. This decryption performance maybe adequate for non-real time systems, particularly if a public key isused to encrypt a secret symmetric-key and send it to the recipientfirst. All subsequent information then can be encrypted using thesymmetric-key, which improves performance, as symmetric-key algorithmsare generally much faster.

In her book, “In Code: A Mathematical Journey”, (ISBN 0-7611-2384-9)Sarah Flannery describes what she calls the “Cayley Purser” public keyalgorithm in Appendix A which requires finding matrices A and C in GL(2,Z_(n)) that are not multiplicatively commutative, i.e.:AC≠CA

The algorithm then requires generating matrix B using:B=C⁻¹A⁻¹C  (A1)

The algorithm further requires generating the matrix G using:G=C^(k)

Where k is a chosen integer greater than 1 or less than −1 so thatmatrix C cannot be trivially found from matrix G. The C matrix is theprivate key. {A,B,G,n} form the public key. The matrix rank is assumedto be 2. In the Postscript of Appendix A [see [6.3], pages 290-292],Flannery describes a security flaw in her algorithm because whencalculating matrix B above, the matrices to the left and right of matrixA in equation (A1) are relatively inverse to each other, so that anylinear multiple of C (modulo n) is also a solution to equation (A1).

In many network applications, client-server models of computerinteractions over networks use context-less servers, where the serverknows nothing about the client, so all context-specific information iskept on client systems. Cookies are an example of client contextinformation, which are kept on client systems instead of web servers.

The original IP (Internet Protocol) packet transmission protocol is asession-less packet transmission protocol used widely on the Internet.Any concept of communications sessions is kept at a higher level, forexample, in applications such as TCP (Transmission Control Protocol).The secure version of IP, called IPSec, is an extremely complexprotocol, designed for all applications requiring use of IP. It istherefore used in a session-less manner, i.e., it is not informed whencommunication sessions begin and end. To minimize the slowness of publickey systems, IPSec frequently uses secret (symmetric) key encryption anddecryption, where the same key is used to both encrypt and decrypt amessage. This in turn requires a secret key exchange, followed bykeeping secret keys at both ends of the secure communications path for aperiod of time that is invisible at the application layer. This secretkey persistence is termed a SA (Security Association). SAs are notinstantiated at the application level, but must occur and be maintainedby IPSec itself, while IPSec is being used in a session-less manner byapplications. This makes maintenance of a security state on amulti-client system such as a web server a very complex task, requiringexpiring and overlapping SAs, and increased use of processor and memoryresources.

To provide context-less servers with public-key encryption, it isdesirable not to keep client-specific private symmetric-keys on theserver. In this case, the slow decryption rate of public keys can be aproblem, even when they are used only to exchange a secret key. Further,the processing requirements for performing simultaneous encryption anddecryption should be reduced, allowing for use in low-powerapplications, such as cell phones, or web-based radio communicationsystems, such as, blue-tooth and wire-less LAN.

Thus, there is a need for a public key system that can perform bothencryption and decryption with relatively fewer calculations, which canresult in a higher encryption/decryption throughput, and/or lower powerconsumption.

BRIEF DESCRIPTION OF THE DRAWINGS

FIG. 1 is a flowchart illustrating one embodiment of generating andpublishing a public key for secure data transmission according to theclaimed subject matter.

FIG. 2 is a flowchart illustrating one embodiment of encrypting data tobe transmitted using the published public key shown in FIG. 1 accordingto the claimed subject matter.

FIG. 3 is a flowchart illustrating one embodiment of decrypting thetransmitted encrypted data shown in FIG. 2 according to the claimedsubject matter.

FIG. 4 is an illustration including a vector diagram of one embodimentof an exchange of secret matrix K between a sender and a receiveraccording to the claimed subject matter.

FIG. 5 is a schematic diagram illustrating an example embodiment of asecure data transmission system according to the claimed subject matter.

FIG. 6 is a block diagram of an exemplary computer system implementingembodiments of the present invention, such as those shown in FIGS. 1-5.

DETAILED DESCRIPTION OF THE INVENTION

In the following detailed description, reference is made to variousspecific embodiments in which the invention may be practiced. Theseembodiments are described with sufficient detail to enable those skilledin the art to practice the invention, and it is to be understood thatother embodiments may be employed, and that structural, logical,electrical, and process changes may be made without departing from theteachings of the invention.

In the foregoing description of the preferred embodiments, variousfeatures of the invention are grouped together in a single embodimentfor the purpose of streamlining the disclosure. This method ofdisclosure is not to be interpreted as reflecting an intention that theclaimed invention requires more features than are expressly recited ineach claim. Rather, as the following claims reflect, inventive aspectslie in less than all features of a single disclosed embodiment. Thus,the following claims are hereby incorporated into the description of thepreferred embodiments, with each claim standing on its own as a separatepreferred embodiment of the invention.

Basic Terms and Notations

Lower case characters in the following description represent integers.Upper case characters represent matrices. { } denotes a comma-separatedset of values. Square matrices form rings under addition andmultiplication because they have the following properties:(A+B)+C=A+(B+C); associative under additionA+B=B+A; commutative under additionA+0=A; the 0 matrix (with rank matching A) forms the additive identityA+(−A)=0; the additive inverse of A is −A. They sum to the 0 matrixA(BC)=(AB)C; associative propertyAI=IA=A; a (multiplicative) identity matrix I exists which commutes withall AA(B+C)=AB+AC; right distributive(A+B)C=AC+BC; left distributive

Matrices form a group under multiplication when, in addition to theproperties of rings above, the following is also true:AA⁻¹=A⁻¹A=I

In other words, for every member of the group A, a unique(multiplicative) inverse matrix exists. Matrices of rank r with integerelements (i.e. elements in Z) are referred to as Mat(r,Z). Matrices withelements in Z, modulo n, are referred to as Mat(r,Z_(n)). Squarematrices of rank r with elements in Z modulo n, and for which an inversematrix exists, are said to belong to the General Linear groupGL(r,Z_(n)).

The (multiplicative) inverse k⁻ of an integer k is calculated, modulo n,(note: the term “reciprocal” is otherwise used when not working in thering of integers modulo n), such that:kk ⁻¹≡1(mod n)

The value k⁻¹ is an integer, and it exists and can be found as long asGCD(n,k)=1, where the well-known GCD( ) function finds the GreatestCommon Divisor of {n,k}. The inverse of k therefore depends on n, verydifferent from the reciprocal of k. The well-known Extended GCD( )algorithm is used to actually find multiplicative inverses. If theextended GCD(n,k) returns a value not equal to 1 (an extremely unlikelychance for large n), then the value is a factor of n. The Extended GCDfunction is also used in matrix inversion, where all elements in theinverse matrix require multiplication by the multiplicative inverse ofthe determinant of the input matrix, modulo n.

The present invention provides techniques for secure data transmissionusing a public key system. An embodiment of a method of the presentinvention is described using FIGS. 1-3. The flowcharts illustrated inFIGS. 1-3 include operations, which are arranged serially in theexemplary embodiment. Many operations in the flowcharts showcalculations of results that depend on other previous results. Anyre-ordering of these calculations in an embodiment which maintains thesedependencies must be viewed as falling under the scope of thisinvention. However, other embodiments of the invention may execute twoor more operations in parallel using multiple processors or a singleprocessor organized as two or more virtual machines or sub-processors.Moreover, still other embodiments may implement the operations as two ormore specific interconnected hardware modules with related control anddata signals communicated between and through the modules, or asportions of an application-specific integrated circuit. Thus, theexemplary process flow is applicable to software, firmware, and hardwareimplementations.

FIG. 1 is a flowchart illustrating one example embodiment of a process100 of generating and publishing a public key for secure datatransmission according to the present invention.

The process begins with operation 110 by finding two unique randomsecret prime numbers p and q. Operation 120 includes generating aninteger modulus n using prime numbers p and q. The modulus n is computedaccording to the equation:n=pq  (1)

Generally, large prime numbers are chosen for p and q to prevent easyfactorization of n to obtain the set of factors {p,q}. Note: if {p,q}are revealed to an attacker, then the roots of integers in Z_(n) can befound rapidly, and at that point the public key is insecure. In theseembodiments, the chosen prime numbers may be discarded after computingmodulus n, or they can be kept to speed up message decryption using thewell-known Chinese Remainder Algorithm.

All matrices described in the present invention belong to the closedLinear Group of matrices GL(r,Z_(n)), unless otherwise stated.

Operation 130 includes generating two matrices {A, C} in GL(r,Z_(n)),such that:AC≠CA  (2)

i.e. matrices {A, C} are a pair of matrices that are notmultiplicatively commutative modulo n. Note that some randomly chosenpairs of matrices are commutative, but this is statistically veryunlikely for large n. For example, a matrix commutes with itself, andtherefore with any matrix which is a power of itself. Non-commutativebehavior is different than integer multiplication. In the later case,the order of multiplication makes no difference to the product obtained.Wherein r is the rank of matrices and Z_(n) denotes elements in integermodulo n.

To ensure that matrices {A, C} are both in GL(r,Z_(n)), both matrixdeterminants should be tested for relative primeness to n (determinedusing a GCD algorithm). Note that 0 is divisible by n, so GCD(0,n)=n andtherefore 0 is not relatively prime to n. The chance of finding eitheri) commutative behavior or ii) common prime factors with n are extremelylow for a large n, so checking for these properties may be omitted.However, applying the checks allow algorithm implementations to bereliably tested for a small n. Smaller rank matrices are preferable(particularly r=2) because encryption and decryption is faster. Higherrank matrices can be used (i.e., r>2), but they result in morecomputation, larger minimum message sizes, and generally no significantimprovement in security because, as is seen later, the best way ofbreaking this type of encryption is to factorize n.

Operation 140 includes generating a matrix B in the group using theequation:B=CAC  (3)

Wherein {A, C} are matrices found in operation 130.

Operation 150 includes generating a matrix G that is multiplicativelycommutative to C, modulo n, i.e. CG=GC. One embodiment of doing thisuses the fact that powers of a matrix commute, so that:G=C^(k)  (4)where k is an integer. To be provably secure, k must be even, andpreferably small (e.g., k=2) for faster key generation. If r=2, then kcannot be an odd number that is small or guessable. In all cases, kcannot belong to the set {−1, 0, 1}.

In another embodiment, G is generated using a linear combination ofpowers of C up to r−1 using:

$\begin{matrix}{G = {\sum\limits_{i = 0}^{r - 1}{u_{i}C^{i}}}} & (5)\end{matrix}$where each of the r values of u_(i) are randomly generated values inZ_(n), and preferably with at least one U_(i)≠0 for i>0 so that G doesnot commute with A. From (5), the number of combinations of u_(i) valuesgives the number of possible G matrices from a given C as n^(r).Generating truly random values in the inclusive range 0 to n−1 ispossible (at a fairly slow rate) on a computer, but hardware basedrandom number generators can give much better number generating rates.As random number generation is also needed for message encryption, thismay be a good alternative method. The generation of random numbers canbe a significant overhead for short messages, so hardware based randomnumber generators may be used to economically generate random sets ofvalues of u and v in equations (5) or sets of v in equation (6).

Operations 160 and 170 include keeping matrix C as the private key, andforming {A,B,G,n,r} as the public key, respectively. In a standardizedalgorithm, a known value of r (e.g., r=2) may be assumed, so r may notneed to be in the public key. Operation 180 includes publishing theformed public key {A,B,G,n,r} for encrypting messages to be transmitted.

FIG. 2 is a flowchart illustrating one embodiment of a process 200 ofencrypting data to be transmitted using the published public key shownin FIG. 1 according to the present invention.

The process begins with operation 210 by obtaining the published key A,B, G, and n, and matrix rank r for encrypting a message to betransmitted. Operation 220 includes generating a first random secretmatrix D that is commutative with the obtained matrix G. In someembodiments, the first random secret matrix D is generated using theequation:

$\begin{matrix}{D = {\sum\limits_{i = 0}^{r - 1}{v_{i}G^{i}}}} & (6)\end{matrix}$where G⁰=I, the identity matrix with the same rank as G, and v_(i) forma set of r secret and independently random integers modulo n, and atleast one v_(i)≠0 for i>0 so that D does not commute with A.

Operation 230 includes generating a second secret key matrix K using thegenerated matrices B and D (later we shall see that K is able to befound easily by the receiver, but not by an eavesdropper). The secondsecret key matrix is computed according to the equation:K=DBD  (7)

Operation 240 includes generating a message matrix E using the receivedpublic key matrix A, and the generated matrix D. The message matrix E isgenerated according to the equation:E=DAD  (8)

Operation 250 includes encrypting a message to be transmitted byapplying all or part of K from (7) as an encryption key in asymmetric-key encryption system. A symmetric-key encryption algorithmuses the same key to encrypt and decrypt a message, so if the messagerecipient can re-obtain K, then decryption is also possible. Examples ofsymmetric key ciphers include DES (Data Encryption Standard), IDEA(International Data Encryption Algorithm), FEAL (Fast Data EnciphermentAlgorithm), RC5, etc.

In one embodiment, a symmetric-key encryption algorithm includespartitioning and packaging an obtained message into a sequence ofunencrypted matrices U_(i). Then each of the unencrypted matrices areencrypted to form a series of corresponding encrypted matrices such thatU_(i)′=KU_(i)K.

Operation 260 includes transmitting the generated message matrix E alongwith the encrypted message. The one or more encrypted matrices must betransmitted in a known index order.

FIG. 3 is a flowchart illustrating one embodiment of a process 300 ofdecrypting the transmitted encrypted message shown in FIG. 2 accordingto the present invention.

The process begins with operation 310 by receiving the transmittedmessage matrix E along with the encrypted message. In some embodiments,operation 310 includes receiving the sequence of encrypted matricesU_(i)′ obtained by using the symmetric-key encryption algorithm.

Operation 320 includes retrieving the second secret matrix K using thereceived message matrix E and the kept private key matrix C. The secondsecret matrix K is calculated according to the equation:K=CEC  (9)

Operation 330 includes decrypting the received encrypted message usingthe retrieved second secret matrix K. In some embodiments, the receivedencrypted message is decrypted by applying all or part of K from (9) asa decryption key in a symmetric-key encryption system. Symmetric-keyencryption uses the same key to encrypt and decrypt a message, so if themessage recipient can re-obtain K, then decryption is also possible. Insome embodiments, the received encrypted message is decrypted byobtaining a matrix Q that is the multiplicative inverse of the retrievedsecond secret matrix K using Q=K⁻¹. Obtained matrix Q is then used todecrypt each of the obtained encrypted matrices U_(i)′ to retrieve thecorresponding decrypted matrices U_(i) by using Q U_(i)′Q. Decryptedmatrices U_(i) are then unpacked and concatenated to obtain thetransmitted message.

Proof that K can be Found by the Receiver

By using (8) to eliminate E in (9), we get:K=CDADC  (10)

From (4) or (5), C and G are multiplicatively commutative, and from (6),G commutes with D. Therefore C commutes with D, so that (10) can berewritten as:K=DCACD  (11)

From (3), we replace CAC with B to obtain:K=DBD  (12)which agrees with (7), proving that (9) correctly obtains K at thereceiver.

Proof of Security of the Public Key

The reason the modulus n is a product of two unknown primes is becauseit can be made extremely difficult to find its prime factors for asufficiently large n, because it is well known that finding factors of nis equivalent to finding roots modulo n. From (4), C may be found from Gif k is known or small. Therefore, the method cannot use a prime numberfor n, as n is then its own prime factorization. The method could usevalues of n with more than two prime factors, but for a given size of n,such values are easier to factorize. Larger values of n require moreaccurate and slower computations. Therefore, for a given computationaleffort, the best security is obtained when two large random primes areused, with roughly comparable sizes. This type of modulus n is also theapproach used in the RSA public key modulus, and some other public keymethods such as Rabin's scheme. As will be seen later, the proof ofsecurity of the new public key algorithm also relies on thecomputational difficulty of finding square roots modulo n.

The Cayley Hamilton Theorem

Let A be an r×r matrix in Mat(r,Z), and p(x) be its characteristicpolynomial. Then the Cayley Hamilton theorem states that p(A)=0. Thecharacteristic polynomial in x of a matrix A is given by:p(x)=Determinant(A−xI)  (14)where I is the identity matrix with the same rank as A (i.e., I=A⁰).Finding the Smallest Sets of Mutually Commutative MatricesIf a_(ij) represents the elements of a matrix A at row i, column j, thenfor a rank r=2 matrix, the characteristic polynomial p(x) is:

$\begin{matrix}{{{Determinant}\begin{bmatrix}{a_{00} - x} & a_{01} \\a_{10} & {a_{11} - x}\end{bmatrix}} = {{\left( {a_{00} - x} \right)\left( {a_{11} - x} \right)} - {a_{01}a_{10}}}} & (15)\end{matrix}$Collecting terms in x on the right, we obtain the characteristicpolynomial as:p(x)=x ² −x(a ₀₀ +a ₁₁)+(a ₀₀ a ₁₁ −a ₀₁ a ₁₀)  (16)Substituting A for x (according to the Cayley-Hamilton theorem) andsetting the result equal to 0 gives:p(A)=A ² −A(a ₀₀ +a ₁₁)+(a ₀₀ a ₁₁ −a ₀₁ a ₁₀)I=0  (17)This result implies that A² can be expressed in the form:A ² =k ₁ A+k ₂ I  (18)i.e. the square of matrix A can be expressed as a linear sum of A andthe identity matrix. This also implies that:A ³ =AA ² =A(k ₁ A+k ₂ I)=k ₁ A ² +k ₂ A  (19)The A² term in the right expression can be eliminated using (18) togive:A ³ =k ₁(k ₁ A+k ₂ I)+k ₂ A=(k ₁ +k ₂)A+k ₁ k ₂ I  (20)

Continuing this procedure, it can be seen that any power of A can bedecomposed into a linear sum of A and I. Similar results can be obtainedfor higher rank matrices, where a matrix A of rank r has acharacteristic polynomial up to degree x^(r) which can be decomposedinto a linear sum of all powers of A from 0 to r−1, e.g., a rank 5matrix of any power can be decomposed into a linear sum of its powersfrom 0 (the identity matrix) to 4. Therefore, for any A of rank r, andinteger power m:

$\begin{matrix}{A^{m} = {\sum\limits_{i = 0}^{r - 1}{u_{i}A^{i}}}} & (21)\end{matrix}$for some set of r values of u_(i). This result shows the equivalencebetween (4) and (5). Any equality in Mat(r, Z) is also true in GL(r,Z_(n)), so if B is a matrix in GL(r, Z_(n)), then the following must betrue:

$\begin{matrix}{B^{m} = {\sum\limits_{i = 0}^{r - 1}{u_{i}B^{i}}}} & (22)\end{matrix}$

This result also shows that any polynomial in B can be decomposed intothis form of sum. All combinations of r values of u_(i) modulo n willthen generate all the members of the commutative set. The number of setmembers is given by:members(n,r)=n ^(r)  (23)members of the commutative set (but not all members of the commutativeset in some cases, as discussed later). This result agrees with thenumber of matrices G that can be generated from all possiblecombinations of u_(i) and a given C in (5). This result is the minimumcommutative set size, and it is the guaranteed minimum number ofmatrices to search for secret matrix C (knowing G—another member of thesame commutative set) by brute force, should this approach be taken tobreak the public key. This is actually a far greater search space than abrute force search for prime factors of n, and is not a feasibleapproach to breaking the public key.

Each set contains all possible multiples of the identity matrix, modulon, so the number of matrices that do not commute with any others outsideof the set is:ExclusiveMembers(n,r)=n ^(r) −n  (24)

For an r×r matrix with elements modulo n, the total number of possiblematrices is:matrices(n,r)=n ^(r×r)  (25)

If the characteristic polynomical of the matrix A is factorizable, thenit can be expressed as a product of a set of lower order polynomials inx. If some product of a subset of these polynomials is zero, then wehave a reduced degree polynomial in A (compared with (21)) that is equalto zero, and the reduced polynomial is no longer uniquely characterizedby the matrix A.

For example, a rank r=3 matrix A will have a cubic characteristicpolynomial with polynomial terms in x. If that polynomial isfactorizable, then it is possible for two of these roots to multiply tozero in the group when A is substituted for x. The existence of reduceddegree polynomials is only possible because of modulus n when working inGL(r, Z_(n)). This matrix A then has a reduced degree polynomial factor.

It is therefore possible for other matrices to have a factorizablecharacteristic polynomial that shares this same reduced degreepolynomial, so the members of these sets will also commute with the setassociated with matrix A. This proves the existence of largercommutative sets than defined in (23) when r>2. However, it is easy toshow that upper triangular or lower triangular matrices do not commutefor any GL(r, Z_(n)), so we know that we cannot chose a group where allmatrices commute. In fact, it can be shown that a minimum bound on thenumber of non-commutative sets is:n^(2(r−1))  (26)

For proof of security of the public key, it is sufficient to show thatthere is a minimum number of members in each commutative set (makingsearches based on a known member of the set impossibly difficult), andthat more than one set exists, allowing large combinations of pairs ofnon-commutative matrices A and C exist in (2).

Given the large number of non-commutative sets, it is also very easy torandomly generate suitable pairs of matrices {A,C} in GL(r,Z_(n)) neededin (2)—in fact, the chance that A and C belong to the same commutativeset is, from (26), 1 in n^(2(r−1)) or less, although this is onlyrelevant when considering the speed of generating public keys, and isnot relevant to public key security.

It is well known that finding the kth roots of G modulo n from (4) aloneis equivalent to factorizing n, which is assumed to be impossiblydifficult for large enough n, even in the simplest non-trivial case whenk=2.

A more effective approach to breaking a key attempts to utilize allknown information about the public key, although in the process of usingthis information, it is then shown that within certain constraints, thekey can be proven to be secure. The following equations apply to any kin (4) with rank r=2 matrices, and we later see that use of either smallodd k or known odd k is insecure. First we define a known M from theknown matrices B and G in the public key:M=BGB⁻¹  (27)

From (3), B contains matrix A as a factor, and from (4), G commutes withC, so from (2) BG≠GB and therefore M≠G. Next we eliminate B from (27)using (3) to get:M=(CAC)G(CAC)⁻¹  (28)

From (4), C and G must commute, so we swap the G with a neighboring C,and expand out the inverse matrices to get:M=CAGCC⁻¹A⁻¹C⁻¹=CAGA⁻¹C⁻¹  (29)

As C commutes with G and not matrix A, then G does not commute withmatrix A. The known matrix N is defined from known public key matrices{A, G} as:N=AGA⁻¹  (30)

As GA≠AG then N≠G. From (29) and (30), we get:M=CNC⁻¹  (31)

N contains matrix A as a factor, so CN≠NC and therefore M≠N. The form of(31) is similar to the CP algorithm public key with the security flaw.The following attack on the new algorithm is based on thistransformation, but unlike the CP algorithm, we later find that theattack applies only for odd, guessable values of k.

The form of (31) allows us to find a linear multiple of C, i.e. uC, butneither of {u, C} are known. However, unlike the CP algorithm, only whenu²≡1 mod n will the congruence (3) be satisfied, as the values of u donot cancel. This is the principal reason for the security of the newalgorithm for rank r=2 matrices.

From (4), G is a power of C, so the result (22) obtained from theCayley-Hamilton theorem allows the definition of G for matrices withrank r=2 to be expressed as a linear combination of the identity matrix(with matching rank r=2) and C as:G=u ₀ I+u ₁ C  (32)

This relationship is implicit using (4) or explicit using (5) in thepublic key generating algorithm, but method (5) generates the equivalentof raising C to a large, unknown power, k, so it is secure. However, itrequires generating truly random numbers, so it is worth proving thesecurity of (4). We now see how it may be possible to find C for rankr=2 matrices under certain circumstances. Note that from (5), higherrank matrices have too many unknowns in u_(i), so they are notvulnerable to this attack. However, higher rank matrices arecomputationally more expensive, so the r=2 case is the most useful tocharacterize. From (32), a linear multiple of C can be obtained in termsof unknown v_(i) values as:v ₁ C=G+v ₀ I  (33)

From (31), multiply both sides by C on the right:MC=CN  (34)

Scaling both sides by v₁:Mv₁C=v₁CN  (35)and substituting for v₁C from (33), we get:M(G+v ₀ I)=(G+v ₀ I)N  (36)

Collecting terms with v_(o)I on the left, and others on the right ({v₀,I} both commute with everything), we get:v ₀ I(M−N)=GN−MG  (37)so thatv ₀ I=(GN−MG)(M−N)⁻¹  (38)

We know that M≠N from (31), so either matrix inversion is possible, orelse n is factorized. From (33), and using (38) to eliminate v₀I, weget:v ₁ C=G+(GN−MG)(M−N)⁻¹  (39)

The right side of (39) consists entirely of known matrices, andtherefore the product v₁C can be found. Multiplying with v₁ iscommutative, so the identityv ₁ ² CACB ⁻¹=(v ₁ C)A(v ₁ C)B ⁻¹  (40)is true, which simplifies on the left using (3), so we can find v₁ ² as:v ₁ ² I=(v ₁ C)A(v ₁ C)B ⁻¹  (41)

From (39), we know v₁C, and we know {A,B} from the public key, so thefactor v₁ ² can be found. To break the key using (41), a square-root ofv₁ ² modulo n has to be found in order to find v₁, and then find C fromthe known v₁C. Obtaining such a square root is known to be equivalent tofactorizing n. Therefore, results (39) and (41) cannot be used as abasis for an attack on their own.

Now we make use of the relationship between C and G in (4):(v ₁ C)^(k) =v ₁ ^(k) C ^(k) =v ₁ ^(k) G  (42)so we can find:v ₁ ^(k)=(v ₁ C)^(k) G ⁻¹  (43)as v₁C is known from (39), and G is known from the public key.

From (41), we know v₁ ², so if k is a finite unknown odd integer, then asearch by repeated division of v₁ ² into v₁ ^(k) from (43) willeventually yield a remaining factor v₁. The value can be rapidlyverified as correct for each search step by squaring the obtained valueand comparing with v₁ ² in (41). If k is a known odd integer, then thenumber of times v₁ ² divided into v₁ ^(k−1) can be immediately found as(k−1)/2. Then v₁ ^(k−1) can be found easily using a power-mod functionbased on successive squarings of v₁. The multiplicative inverse of v₁^(k−1) is then obtained (modulo n) and multiplied by v₁ ^(k) from (43)to obtain v₁. Once v₁ is found, then C can be found from (39), and thepublic key is broken. If k is a large, unknown odd integer modulo n,then the key is secure against this particular attack, but its securityhas not been proved.

In the Cayley-Purser public key algorithm, generating the public key bycalculating C to the power k modulo n in (4) is a relatively slowoperation, particularly compared to multiplication for large k, so forefficient implementations with rank 2 matrices, k should be relativelysmall. For the public key to be secure in this case, k must be an even,non-zero integer. k=2 is the fastest value for calculating (4) that isalso secure from this attack. For higher rank matrices, the attack doesnot apply, but the same value of k=2 also works well.

Having proven that G can be generated using the CP public key algorithm,it has been shown that the code is not secure for rank 2 matrices andfor odd k that is known or guessable. We now prove the public key issecure for known even k. We do this by creating a public key accordingto (1) . . . (3), and by generating a second public key derived from thefirst one. We then show that finding C for both keys is equivalent tofinding the √{square root over (d)} modulo n for a chosen d. In thederived public key, we assume the existence of a new secret matrix C′that is related to C by:C′=√{square root over (d)}C  (44)and that both C′ and √{square root over (d)} are unknown. Also, {A,n}are shared between the C and the derived C′ public keys. Substituting C′for C in (3), we get B′ of the derived key as:B′=C′AC′  (45)and eliminating C′ using (44), and collecting √{square root over (d)}terms, we get:B′=dCAC=dB  (46)

Note that B′ is calculated from {B,d,n}, and not from √{square root over(d)} or C′, which are unknown when the key is generated.

By also substituting C′ for C in (4), and using the fact that k is evenso k/2 is an integer, we get:G′=(C′ ²)^(k/2)  (47)Then we eliminate C′ using (44) to obtain:G′=(dC ²)^(k/2)  (48)By regrouping powers:G′=d ^(k/2) C ^(k) =d ^(k/2) G  (49)

Note that G′ is calculated from {G,k,d,n} only (and not from √{squareroot over (d)} and C′ which are unknown).

The following steps now complete the proof:

-   -   Step (a): A first public key is generated according to (1) . . .        (4), and with a large composite n whose factors are unknown, so        the public key {A,B,G,n} and the exponent k and private key        matrix C are all known.    -   Step (b): A second public key (related to the public key        generated at Step (a)) is generated. First, a value d is chosen        for which the value of √{square root over (d)} is unknown        modulo n. The new public key consists of {A,B′,G′,n} where B′ is        calculated from B using (46), and G′ is calculated from G        using (49) and the known k. The values {A,n} are the same as in        the key generated in Step (a). Note that neither the matrix C′        nor √{square root over (d)} are used in generating either public        key.    -   Step (c): The matrix C′ is “found” from all available        information—{A,B,B′,C,G,G′,d,k,n} using a hypothetical        polynomial-time algorithm. Note: if C′ cannot be found from        these, then it certainly cannot be found from the derived public        key {A,B′,G′,n}, which is a subset of the above set.    -   Step (d): The value of √{square root over (d)} can now be found        from (44) using C generated in Step (a), and C′ found in        Step (c) using:        √{square root over (d)}I=C′C ⁻¹  (50)    -   Step (e): The value of √{square root over (d)} was not involved        in generating B′ or G′, so finding √{square root over (d)}        modulo n by finding C′ is equivalent to finding √{square root        over (d)} modulo n for a chosen d. If C′ can be found in        polynomial time, then this is computationally equivalent to        factorization in polynomial time, which is assumed to be        impossible. Therefore, as all other steps are simple to perform,        it is not possible to find C′ in Step (c) in polynomial time,        and no hypothetical algorithm for rapidly breaking the public        key exists.    -   QED

If the new public key could be broken, and factorization of n could beachieved from Step (c), then the technique could be used to break RSApublic keys, and Rabin's public key scheme, amongst others.

A Visual Analogy to Obtaining K at the Sender and Receiver

FIG. 4 is a vector diagram 400 that illustrates an analogy of how secretmatrix K is exchanged between a sender and a receiver according to theclaimed subject matter. This diagram shows how, starting from thetop-left at matrix A, both the sender and receiver can arrive at thesame matrix K on the bottom right, by following two different pathsaround a parallelogram. The lengths and directions of each side of theparallelogram represent operators applied to one matrix to obtainanother.

FIG. 4 shows the C( )C operator as two vectors representing the shortersides 415 and 435 of the parallelogram, and the D( )D operator as twovectors representing the longer sides 422 and 424. The important thingto notice is that the vectors on opposite sides of the parallelogram arethe same direction and length, although their depicted directions andlengths are for illustrative purposes only. In actuality, the operatorsare multiplicative rather than additive, so FIG. 4 represents a kind of‘log’ of the operators, allowing vector length addition to be used.Also, the present invention uses matrices in the group GL(r, Z_(n)),which cannot actually be rendered onto a two-dimensional surface.

FIG. 4 illustrates a far left portion 410, a middle portion 420, and afar right portion 430. The far left portion 410 of the vector diagram400 illustrates using a vector 415 to represent operator C( )C as afunction applied to A to obtain B=C(A)C as in (3). The length anddirection of the vector represents the scaling affect of the C( )Coperator, and is uniquely determined by secret matrix C.

The middle portion 420 of the vector diagram 400 assumes the sender hasobtained the published public key transmitted by the receiver, and isusing it to generate a first random secret matrix D that is commutativewith the kept matrix C and generated matrix G. The sender does not knowthe direction of the operator vector C( )C, but does know the two pointsin the parallelogram at {A,B} from the published public key. The messagesender can then apply the D( )D operator to the matrices {A,B} to obtainE=D(A)D from (8), and K=D(B)D from (7), represented by scaling androtating vectors 422 and 424 in FIG. 4. Note that the sender obtains Kby traversing first the C( )C path from the top left at matrix A to B,and then the D( )D path from B to K. The sender then encrypts messagesto be transmitted using the generated second secret matrix K andtransmits the encrypted messages along with the message matrix E.

The far right portion 430 of the vector diagram 400 illustrates thereceiver receiving the transmitted encrypted messages along with themessage matrix E from the sender. Vector diagram 400 further illustratesusing a vector 435 how, from the sender in the middle portion 420, thereceiver obtains the message matrix E via the D( )D path across the topof the parallelogram, from the starting point at matrix A. As thereceiver is also the public key generator, the C matrix is known. The C()C operator can then be applied to the E matrix to obtain the final pathdown to K. The Receiver's route to K from A is therefore via the D( )Dto E, and then the C( )C operator. This route around the parallelogramis different from the route that the sender took from A to K describedearlier. This is equivalent to saying that the C( )C and D( )D operatorscommute with each other—in other words, they can be applied in eitherorder from A to K. Also note that the receiver manages to obtain Kwithout knowing D, and similarly, the sender gets from A to K withoutknowing C.

FIG. 5 illustrates one embodiment of a system 500 used for secure datatransmission according to the present invention. FIG. 5 includes anexample of a system 500 having a sender 510 coupled to a network 520. Inaddition, FIG. 5 includes a receiver computing platform 530 coupled tonetwork 520. Further, FIG. 5 includes a receiver 540 coupled to bothnetwork 520 and receiver computing platform 530. In some embodiments,sender 510, receiver computing platform 530, and receiver 540 arecoupled to network 520 through a transmission medium 550. Sender 510includes an encoder 515. Receiver computing platform 530 includes memory532 and a processor 534. Receiver 540 includes a decoder 545.Transmission media 550 may include, for example, fiber optic cable,category 5 (CAT-5) networking cabling, or wireless media such aswireless local area network (LAN).

In operation, processor 534 generates two matrices A and C of rank r andwith each element in the integers modulo n such that AC does not equalCA. In these embodiments, modulo n is obtained as a product of twounique randomly chosen secret prime number p and q. Also, in theseembodiments, r is the rank of the matrices A and C. Processor 534 thengenerates matrix B by using the generated matrices A and C such thatB=CAC. Process 534 further generates matrix G such that the generatedmatrix G is in the same multiplicatively commutative subgroups as matrixC. Processor 534 then publishes the generated matrices A, B, G, andmodulo n and matrix rank r as the public key and retains the generatedmatrix C as the private key. The generation of matrices A, B, C, G, andmodulo n and matrix rank r are explained in more detail with referenceto FIG. 1. In some embodiments, memory 534 stores the generated publicand private keys.

Encoder 515 desiring to transmit a secure message obtains the publishedpubic key including matrices A, B, G, and modulo n and matrix rank rthrough the network 550 to encrypt the message to be transmitted.Encoder 515 then generates a first random secret matrix D that ismultiplicatively commutative with the obtained matrix G. Encoder 515then generates a second secret matrix K using the generated matrices Band D. Encoder 515 then obtains a message to be transmitted securely andencrypts the obtained message using a symmetric-key algorithm by usingthe generated secret matrix K as the key. Encoder 515 then transmits thegenerated message matrix E along with the encrypted message. Thegeneration of first and second secret matrices D and K are explained inmore detail with reference to FIG. 2.

Decoder 545 desiring to receive the encrypted message receives thetransmitted message matrix E and the encrypted message through thenetwork 550. Decoder 545 then retrieves the second secret matrix K usingthe received message matrix E and the privately kept matrix C. Decoder545 then decrypts the received encrypted message using the retrievedsecond secret matrix K with a symmetric-key decryption algorithm, toobtain the transmitted message. The process of using symmetric-keyencryption and decryption to encrypt and decrypt messages, respectively,is explained in more detail with reference to FIGS. 2 and 3.

FIG. 6 is a block diagram of a system according to one embodiment of thepresent invention. Computer system 600 contains a processor 610 and amemory system 602 housed in a computer unit 605. Computer system 600 isbut one example of an electronic system containing another electronicsystem, e.g., memory system 602, as a subcomponent. The user interfacecomponents include a keyboard 620, a pointing device 630, a monitor 640,a printer 650, and a bulk storage device 660. It will be appreciatedthat other components are often associated with computer system 600 suchas modems, device driver cards, additional storage devices, etc. It willfurther be appreciated that processor 610 and memory system 602 ofcomputer system 600 can be incorporated on a single integrated circuit.Such single-package processing units reduce the communication timebetween the processor and the memory circuit. Any of these components ofthe system may contain a memory device that stores instructions that canbe executed by a processor to perform the secure data transmission ofthe present invention.

The above description illustrates preferred embodiments, which achievethe features and advantages of the present invention. It is not intendedthat the present invention be limited to the illustrated embodiments.Modifications and substitutions to specific process conditions andstructures can be made without departing from the spirit and scope ofthe present invention. Accordingly, the invention is not to beconsidered as being limited by the foregoing description and drawings,but is only limited by the scope of the appended claims.

1. A method, comprising: obtaining a message; retrieving a publishedpublic key including matrices A, B, G, and n and matrix rank r forencrypting the message, wherein matrix B is generated by using generatedmatrices A and C such that B≡CAC; generating a first random secretmatrix D that is multiplicatively commutative with the matrix G;generating a second secret matrix K such that K=DBD; generating amessage matrix E such that E=DAD; encrypting the message to provide anencrypted message using a symmetric-key algorithm by using the secondsecret matrix K as a key; and transmitting the encrypted message and themessage matrix E.
 2. The method of claim 1, wherein the encryptedmessage and the message matrix E are sent together to at least onerecipient.
 3. The method of claim 1, wherein encrypting the messageusing a symmetric-key algorithm further comprises: partitioning andpacking the message into a sequence of unencrypted matrices U_(i); andencrypting each of the unencrypted matrices to form a sequence ofcorresponding encrypted matrices such that U_(i)′=KU_(i)K.
 4. The methodof claim 1, wherein generating the first random secret matrix D that ismultiplicatively commutative with the matrix G, comprises: generatingthe first random secret matrix D using the equation:$D = {\sum\limits_{i = 0}^{r - 1}{v_{i}G^{i}}}$ wherein G⁰=I, theidentity matrix having a rank equal to a rank of the matrix G, and v_(i)form a set of r secret and an independently random integers modulo n,and at least one v_(i)≠0 for i>0 so that D does not commute with A.
 5. Acomputer-readable medium having computer-executable instructions, thecomputer-executable instructions comprising: obtaining a message to besent securely; retrieving, from a data store, the published public key,the published public key including matrices A, B, G, and n forencrypting a message to be transmitted, wherein each of the retrievedmatrices have a predetermined matrix rank of r, wherein matrix B isgenerated by using generated matrices A and C such that B≡CAC;generating a first random secret matrix D that is multiplicativelycommutative with the retrieved matrix G; generating a second secretmatrix K such that K=DBD; generating a message matrix E such that E=DAD;encrypting the message using a symmetric-key algorithm by using thegenerated second secret matrix K as the key; and transmitting thegenerated message matrix E along with the encrypted message.
 6. Thecomputer-readable medium of claim 5, wherein encrypting the messageusing a symmetric-key algorithm further comprises: partitioning andpacking the message into a sequence of unencrypted matrices U_(i); andencrypting each of the unencrypted matrices such that U_(i)′=KU_(i)K. 7.The computer-readable medium of claim 6, wherein the predeterminedmatrix rank r is
 2. 8. The computer-readable medium of claim 6, furthercomprising: receiving the transmitted message matrix E; retrieving thesecond secret matrix K using the received message matrix E and a keptprivate key matrix C; and decrypting the encrypted message using theretrieved second secret matrix K to obtain the message.
 9. Acryptographic communication system, comprising: a receiver computingplatform comprising: a processor to generate two matrices A and C ofrank r with each matrix element in the integers modulus n such that ACdoes not equal to CA, wherein n is obtained by using two unique randomlychosen secret prime numbers p and q, wherein r is a matrix rank of thetwo matrices A and C, wherein the processor generates a matrix B byusing the generated matrices A and C such that B≡CAC, wherein theprocessor further generates a matrix G such that the generated matrix Gis in the same multiplicatively commutative subgroup as C, and whereinthe processor publishes the generated matrices A, B, G, and n and matrixrank r as the public key, and retains the generated matrix C as theprivate key; a sender coupled to the receiver computing platform througha network comprising: an encoder to obtain the published matrices A, B,G, and n and matrix rank r for encrypting a message to be transmitted,wherein the encoder generates a first random secret matrix D that ismultiplicatively commutative with the obtained matrix G, wherein theencoder generates a second secret matrix K=DBD using the generatedmatrices B and D, wherein the encoder generates a message matrix E=DADusing the generated matrices A and D, wherein the encoder obtains amessage to be sent securely and encrypts the obtained message using asymmetric-key algorithm by using the generated second secret matrix K asthe private key, and wherein the encoder transmits the generated messagematrix E along with the encrypted message; and a receiver coupled to thesender through the network and further coupled to the receiver computingplatform, comprising: a decoder to receive the transmitted messagematrix E and the encrypted message, wherein the decoder retrieves thesecond secret matrix K and wherein the decoder decrypts the encryptedmessage uses the retrieved second secret matrix K with a symmetric-keydecryption algorithm to obtain the transmitted message.
 10. The systemof claim 9, wherein the matrices A, B, C, and G are mathematicallydescribed as belonging to a general linear group:GL(r,Z_(n)) wherein r is the matrix rank, and each matrix element is inthe integers modulus n, and where an inverse matrix exists for eachmatrix in the group, and for which the group is closed under operationsof matrix multiplication.
 11. The system of claim 9, wherein the matrixrank r is greater than or equal to
 2. 12. The system of claim 9, whereinthe processor generates the matrix G using the equation: G≠C^(k) whereink is a non-zero, even integer.